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From Classical Extensional Higher-Order Tableau to Intuitionistic Intentional Natural Deduction

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From Classical Extensional Higher-Order Tableau to Intuitionistic Intentional Natural Deduction From Classical Extensional Higher-Order Tableau to Intuitionistic Intentional Natural Deduction Chad E. Brown1 and Christine Rizkallah2 1 Saarland University, Saarbrücken, Germany [email protected] 2 Max-Planck-Institut für Informatik, Saarbrücken, Germany [email protected] Abstract We define a translation that maps higher-order formulas provable in a classical ex- tensional setting to semantically equivalent formulas provable in an intuitionistic inten- sional setting. For the classical extensional higher-order proof system we define a Henkin- complete tableau calculus. For the intuitionistic intensional higher-order proof system we give a natural deduction calculus. We prove that tableau provability of a formula im- plies provability of a translated formula in the natural deduction calculus. Implicit in this proof is a method for translating classical extensional tableau refutations into intuitionistic intensional natural deduction proofs. 1 Introduction We describe a way of translating a simply-typed higher-order formula s into a semantically equivalent formula s0 such that if s is provable in a classical extensional higher-order tableau calculus, then s0 is provable in an intuitionistic intentional higher-order natural deduction cal- culus. A potential application of such a translation is mapping refutations found by a classical extensional tableau-based automated theorem prover like Satallax [4] to corresponding proof terms in an intuitionistic intensional natural deduction based system like Coq [15, 3]. The problem of translating classical logic into intuitionistic logic has a long history. A result by Glivenko [11] states that if s is a propositional tautology, then ::s is intuitionistically provable. This result does not hold for first-order formulas. Kuroda [14] showed a similar result if one translates by double negating the formula and adding double negations to the bodies of universal quantifiers. We recently proved that the Kuroda translation generalizes to give a translation from classical intentional higher-order logic to intuitionistic intentional higher order logic [5]. The generalized Kuroda translation does not suffice in the presence of functional extensionality, for an example see [5]. The work of Gandy [9] provides a method to translate from a higher-order logic with extensionality into a higher-order logic without extensionality. The translation we give in this paper uses ideas similar to those of Kuroda and Gandy to deal with both issues at once. The Gandy translation translates equality with the help of a binary relation and a predicate that are defined by mutual recursion. Our translation is simpler in that it translates equality with the help of a single binary relation that is defined inductively on types. Many logical systems of quite different character are commonly referred to as higher-order logics. Some forms of higher-order logic allow classical reasoning while others only allow in- tuitionistic reasoning. For example, formulas such as 8p:p _:p and 8p:::p ! p are provable if and only if the higher-order logic is classical. Likewise, some forms of higher-order logic allow extensional reasoning while others do not. Formulas such as 8fg:(8x:fx = gx) ! f = g J.C. Blanchette, J. Urban (eds.), PxTP 2013 (EPiC Series, vol. 14), pp. 27–42 27 From Classical Extensional to Intuitionistic Intentional C. E. Brown, C. Rizkallah and 8pq:(p ! q) ^ (q ! p) ! p = q are provable if the higher-order logic is extensional and are not provable otherwise. The formula (λx.::x) = (λx.x) is only provable if the logic is both classical and extensional. The extensionality properties can be factored into propositional extensionality (i.e., boolean extensionality) and functional extensionality (which can itself be factored into extensionality properties η and ξ) [1]. In this paper we consider a tableau system with full extensionality and a natural deduction system with no extensionality. Many automated and interactive theorem provers (e.g., Isabelle/HOL [16], LEO-II [2] and Satallax [4]) are based on classical extensional versions of higher-order logic. Automated and interactive theorem provers are used to create proofs, while the task of a proof checker is checking the correctness of proofs. In order to check a proof, the proof must be represented as an explicit object. A well-known way of representing proofs is via the Curry-Howard-de Bruijn correspondence [13, 8]: a natural deduction proof of P can be encoded as a λ-term of type P . It is easy to write a natural deduction system for higher-order logic for which one can obtain proof terms in an obvious way. However, this creates a natural deduction system which is intuitionistic (not classical) and intentional (not extensional). One way to resolve this mismatch is to add classical extensional axioms to the natural deduction system. Currently, Satallax produces Coq proof terms under the assumption that one has added such axioms to Coq [20]. Using the ideas in this paper, one can avoid assuming such axioms in Coq and still produce Coq proof terms from Satallax refutations (unless Satallax makes use of a choice operator). The basic definitions and lemmas used in the translation have all been formalized and proven in Coq, as described in [18, 19]. However, there is currently no implementation of the translation as a whole. In particular, Satallax has not yet been extended to produce Coq proof terms using these definitions and lemmas. Since the translation changes the formula, a Coq user could not typically call a classical extensional theorem prover like Satallax to request a proof term for an arbitrary formula in Coq. On the other hand, one could use the translation to automatically create a Coq development from a development in a classical extensional simple type theory. The rest of the paper is organized as follows. In Section 2 we give a brief overview of simply typed λ-calculus and in Section 3 we introduce a tableau calculus T . We then present the targeted natural deduction calculus N in Section 4. In Section 5 we present our translation. We prove that it maps T -refutable branches to N -refutable contexts in Section 6. We conclude in Section 7 and provide suggestions for future work. More details about the translation in this paper and other translations can be found in the Master’s thesis of one of the authors [18]. 2 Simply Typed Lambda Calculus 2.1 Syntax We describe simply typed λ-calculus in the style of Church [7]. The set of types T is given inductively: o and ι are types and στ is a type whenever σ and τ are types. The types o (the type of propositions) and ι (the type of individuals) are called base types. Types of the form στ are called function types. Often the function type στ is written as σ ! τ, but we use the shorter notation since there are only base types and function types. We use σ and τ to range over types. For each type σ, let Nσ be an infinite set of names of type σ. Some of the names are logical constants: 28 From Classical Extensional to Intuitionistic Intentional C. E. Brown, C. Rizkallah •> and ? are (distinct) logical constants in No. •^ , _, and ! are (distinct) logical constants in Nooo. σ • For each σ, = is a logical constant in Nσσo. σ σ • For each σ, 8 and 9 are (distinct) logical constants in N(σo)o. The remaining names are variables . Let Cσ be the set of logical constants of type σ and Vσ be S S S the (infinite) set of all variables of type σ. Let N = σ Nσ, C = σ Cσ and V = σ Vσ. 0 C0 For any C ⊆ C we define a family of sets of terms Λσ for each type σ by induction. C0 • For every x 2 Vσ, x 2 Λσ . 0 C0 • For every c 2 Cσ \C , c 2 Λσ . C0 C0 • For every x 2 Vσ and s 2 Λτ , (λx.s) 2 Λστ . C0 C0 C0 • For every s 2 Λστ and t 2 Λσ , (st) 2 Λτ . C0 We defined Λσ relative to a set of logical constants. There are two particular sets of logical constants of interest in this paper: the full set C of logical constants and the set C− := f!; 8σjσ is a typeg: C − C− C0 To simplify notation, we define Λσ to be Λσ and Λσ to be Λσ . Note that Λσ ⊆ Λσ for any 0 S C ⊆ C. An element of Λσ is called a term of type σ.A term is an element of σ Λσ. Terms of the form (λx.s) are called λ-abstractions. Terms of the form (st) are called applications.A formula is a term of type o. We write :s for ((! s)?). We write stu for (st)u, except that :st means :(st). We use the infix notations s ^ t, s _ t, s ! t and s =σ t as shorthand for ^st, _st, ! st and =σ st, respectively. We also write s 6=σ t for :(s =σ t). Using infix notation note that :s is the same term as s !?. We sometimes write 8x : σ:s and 9x : σ:s for 8σ(λx.s) and 9σ(λx.s), respectively. We may also omit the type entirely from a quantified formula, equation or disequation and write 8x:s, 9x:s, s = t or s 6= t when the types are clear from the context. If s is a term, x 2 Nσ and t is a term of type σ, then s[x := t] is defined to be the result of substituting t for x in s via a capture-avoiding substitution. A simultaneous substitution θ substitutes several variables simultaneously. We use the notation θ; [x := t] to mean the simultaneous substitution that agrees with θ on all variables except (possibly) x which is mapped to t.
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